Brief itemsreleased by Linus on April 5. It contains a fair number of fixes. Says Linus: "We should be getting close to a 2.6.21 release, so please update any regression reports you've done."
A few dozen patches have been merged into the mainline git repository since -rc6 was released. Your editor guesses that one more -rc will be needed before 2.6.21 is done.
The current -mm tree is 2.6.21-rc6-mm1. Recent changes to -mm include a number of tweaks for Sony laptops, an enlarged set of paravirt_ops hooks, a new set of /proc files for learning about process memory, a rework of the NFS file locking code, and the signalfd() patches. Andrew notes that -mm is now a "rather large" 25MB patch against the mainline.
The current stable 2.6 kernel is 126.96.36.199, released on April 6; 188.8.131.52 had been released moments earlier. The two patches contain a fair number of fixes, including one for a remotely exploitable crash in the Appletalk code.
For older kernels: 184.108.40.206-rc1 was released on April 11 with about a dozen fixes.
Kernel development news
Internally, the core building block for the device model is the kobject. Objects represented in sysfs - devices, for example - each contain a kobject which, among other things, is the focal point for sysfs access. The kobject also contains a reference count for the containing object which is used to manage its lifecycle. A given kobject and its containing data structure can be deleted when the reference count goes to zero - and not before. Reference counting works, but it can lead to surprises.
As an example, consider a USB device - a mouse, say. When this device is plugged into the system, a suitable device structure (containing a kobject) is created and registered with the kernel. When the mouse is unplugged, that structure is released. But imagine what happens if a user-space process opens a sysfs file associated with the mouse device while it is present, and keeps that file open long after the physical device goes away. The kernel must be able to handle operations on that open sysfs file, even though the driver thinks that the device it represents is long gone. The reference counting in the kobject makes this work - most of the time. The potential for confusion is high, though, especially with drivers which have not been written with this sort of lifecycle management in mind.
Back at the end of March, Tejun Heo posted a discussion of device model lifecycle issues which points out this problem and a few others. His argument is that the need to manage objects with different lifecycles makes programming with the device model hard - something developers have known for some time. Even the core device model maintainers will admit that it's easy to get things wrong.
More recently, Tejun has followed up with a patch set which attempts to simplify the situation. There is a great deal of cleanup work in these patches, and one small API change, but the core change is this: it enables a clean separation of the lifecycles of sysfs objects and the underlying data structures they represent. As a result, it is no longer necessary for code outside of sysfs to be concerned about the fact its data structures may have a shorter life than the sysfs objects representing those structures.
A sysfs directory (which represents a kobject) is represented within the kernel by struct sysfs_dirent. In current kernels, if the sysfs_dirent structure exists, its underlying kobject is expected to exist as well. It is not possible for the kobject to go away as long as the sysfs_dirent structure exists; that means that the structure containing the kobject must continue to exist as long as any references to the sysfs files exist. Tejun's patch works by eliminating that requirement.
In the modified sysfs, each sysfs_dirent contains a new counter called s_active. This counter tracks the number of active references to the object; these references are the ones which involve the associated kobject at the current moment. A user-space process which is holding a sysfs file open will not increase the s_active count until it performs an actual operation on that file, and the reference remains only for as long as it takes to complete the operation. Since most sysfs operations are quite fast, active references will not normally be held for long.
The active count, as it happens, is maintained with an rwsem - a reader/writer semaphore. Active references are tracked as readers, so there can be any number of them outstanding at a given time. The code to obtain an active reference works with a call to down_read_trylock(), meaning that it will take a "lock" (a reference) if one is available, but it will not block if the operation fails. All of the relevant sysfs operations have been changed to obtain active references before referencing the kobject - and they make sure that the reference was granted. If an attempt to obtain an active reference fails, sysfs fails the higher-level operation with -ENODEV.
The only way down_read_trylock() will fail is if another thread holds a writer lock on the semaphore - or is in the process of waiting for the readers to get out of the way so it can get that lock. Should something happen which causes the underlying kobject to go away, the cleanup code will call down_write() on the s_active rwsem in the sysfs_dirent entry, thus taking a writer lock. This call will cause any future attempts to obtain an active reference to fail; it will also block until all currently-existing active references are released.
The end result of all this is that, once the final kobject_put() call has completed for a given kobject, there will be no further attempts to access that kobject from sysfs. The kobject (and its containing data structure) can be safely deleted, and the driver need worry no more about it.
As an added bonus, there is no longer any need to increase module reference counts when sysfs attributes are being accessed. A driver which is being unloaded will release all of its devices, meaning that sysfs will no longer make any calls into the driver module anyway; the module reference count becomes superfluous. So Tejun's patch removes the owner field from attribute structures - a change which ripples through a significant amount of driver code.
There have been some comments on how the patches are implemented, but no disagreement with the ultimate goal; these changes could go in as soon as 2.6.22. Tejun would appear to have more improvements in mind, but, even with no further changes, the current patches go a long way toward making sysfs safer and easier to work with.
To begin, they found out that they could not even boot a kernel with the default configuration. Linux systems normally have a limit of 32768 active processes at any given time. Anybody who has run "ps" will have noted that kernel threads are taking up an increasing number of those slots; your editor's single-processor desktop is running 39 of them. In fact, there are now enough kernel threads on a typical system that they will fill that entire space - and more - on a 4096-CPU machine. This problem is relatively easy to take care of by raising the limit on the number of processes. But it gets more interesting from there.
The init process is the parent of last resort for every other process on the system, including kernel threads. So, on a big system, init has a lot of child processes. These children live on a big linked list; that list must be searched by various functions, including the variants of wait(). If the process being searched for is toward the end of the list, that search can take a long time. Since (1) most kernel threads are long-lived, and (2) new processes are put at the end of the list, chances are that a search will, indeed, be looking for a process at the end.
Then, for the ultimate in fun, load a module into the kernel. The module loading process calls stop_machine_run() when the new module is being linked in; this function creates a high-priority kernel thread for each processor on the system. That thread will grab its assigned CPU and simply sit there until told to exit; while all CPUs are locked up in this way the linking process can be performed. Calling a function like stop_machine_run() is a somewhat antisocial act in the best of times. But, in the 4096-processor system, stop_machine_run() will create 4096 threads, each of which goes on the end of init's child list, and each of which must be searched for when the time comes to clean it up. The result is a system which simply stops for an extended period of time.
One could argue that people with systems that large simply should not load modules, but there is a possibility of pushback from the user community. So other solutions need to be found. Robin Holt's problem report included a simple patch which moves exiting processes to the beginning of the child list. This change solves the immediate problem by making searches for those children find them without having to iterate through all of the long-lived processes which are not going anywhere.
Linus had a couple of alternatives. One was to create a separate list for zombie processes, eliminating that search altogether. Another was to stop making kernel threads be children of the init process since they have little to do with user space in any case. But some developers feel that the real solution might be to start cutting back on the number of kernel threads.
The biggest culprit for kernel thread creation will certainly be workqueues, which, by default, create one thread for every CPU on the system. There are situations which can benefit from multiple threads and CPU locality, but there are undoubtedly many places where all of those threads are not needed. Cleaning them up would help to solve some of the scalability issues; as an added bonus it would remove some of the clutter from ps listings.
In many cases, a workqueue may not be necessary at all. Instead, kernel subsystems could just use the "generic" keventd workqueue (which runs as the events/n threads). There are some issues with using keventd, including indeterminate latency and a small possibility of deadlocks, but, for many situations, it may work well enough.
In other cases, using a thread makes sense. Tasks involving long delays are one example; running a function with multi-second delays in keventd is considered impolite. Work requiring complicated context also benefits from its own thread. But, in a number of cases, those threads need not be created until there is actually some work to be done. A quick ps run on most systems will show threads related to error handling, asynchronous I/O, bluetooth, and more. In the current scheme, they are created at boot (or module load) time and many of them may never do any real work before the system shuts down. Thread creation is cheap, so many of these threads could be created on demand when they are needed.
There are probably some real improvements to be made in this area; all that's needed is somebody with the time and motivation to do the work. In the mean time, those of you with 4096-way systems may need to apply a patch or two.
Christoph Lameter is one of those people for whom the slab allocator does not work quite so well. Over time, he has come up with a list of complaints that is getting impressively long. The slab allocator maintains a number of queues of objects; these queues can make allocation fast but they also add quite a bit of complexity. Beyond that, the storage overhead tends to grow with the size of the system:
Beyond that, each slab (a group of one or more continuous pages from which objects are allocated) contains a chunk of metadata at the beginning which makes alignment of objects harder. The code for cleaning up caches when memory gets tight adds another level of complexity. And so on.
Christoph's response is the SLUB allocator, a drop-in replacement for the slab code. SLUB promises better performance and scalability by dropping most of the queues and related overhead and simplifying the slab structure in general, while retaining the current slab allocator interface.
In the SLUB allocator, a slab is simply a group of one or more pages neatly packed with objects of a given size. There is no metadata within the slab itself, with the exception that free objects are formed into a simple linked list. When an allocation request is made, the first free object is located, removed from the list, and returned to the caller.
Given the lack of per-slab metadata, one might well wonder just how that first free object is found. The answer is that the SLUB allocator stuffs the relevant information into the system memory map - the page structures associated with the pages which make up the slab. Making struct page larger is frowned upon in a big way, so the SLUB allocator makes this complicated structure even more so with the addition of another union. The end result is that struct page gets three new fields which only have meaning when the associated page is part of a slab:
void *freelist; short unsigned int inuse; short unsigned int offset;
For slab use, freelist points to the first free object within a slab, inuse is the number of objects which have been allocated from the slab, and offset tells the allocator where to find the pointer to the next free object. The SLUB allocator can use RCU to free objects, but, to do so, it must be able to put the "next object" pointer outside of the object itself; the offset pointer is the allocator's way of tracking where that pointer was put.
When a slab is first created by the allocator, it has no objects allocated from it. Once an object has been allocated, it becomes a "partial" slab which is stored on a list in the kmem_cache structure. Since this is a patch aimed at scalability, there is, in fact, one "partial" list for each NUMA node on the system. The allocator tries to keep allocations node-local, but it will reach across nodes before filling the system with partial slabs.
There is also a per-CPU array of active slabs, intended to prevent cache line bouncing even within a NUMA node. There is a special thread which runs (via a workqueue) which monitors the usage of per-CPU slabs; if a per-CPU slab is not being used, it gets put back onto the partial list for use by other processors.
If all objects within a slab are allocated, the allocator simply forgets about the slab altogether. Once an object in a full slab is freed, the allocator can relocate the containing slab via the system memory map and put it back onto the appropriate partial list. If all of the objects within a given slab (as tracked by the inuse counter) are freed, the entire slab is given back to the page allocator for reuse.
One interesting feature of the SLUB allocator is that it can combine slabs with similar object sizes and parameters. The result is fewer slab caches in the system (a 50% reduction is claimed), better locality of slab allocations, and less fragmentation of slab memory. The patch does note:
Causing bugs to stand out is generally considered to be a good thing, but wider use of the SLUB allocator could lead to some quirky behavior until those new bugs are stamped out.
Wider use may be in the cards: the SLUB allocator is in the -mm tree now and could hit the mainline as soon as 2.6.22. The simplified code is attractive, as is the claimed 5-10% performance increase. If merged, SLUB is likely to coexist with the current slab allocator (and the SLOB allocator intended for small systems) for some time. In the longer term, the current slab code may be approaching the end of its life.
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